Managing speculative memory access requests in the presence of transactional storage accesses

ABSTRACT

In at least some embodiments, a cache memory of a data processing system receives a speculative memory access request including a target address of data speculatively requested for a processor core. In response to receipt of the speculative memory access request, transactional memory logic determines whether or not the target address of the speculative memory access request hits a store footprint of a memory transaction. In response to determining that the target address of the speculative memory access request hits a store footprint of a memory transaction, the transactional memory logic causes the cache memory to reject servicing the speculative memory access request.

PRIORITY CLAIM

This application is a continuation of U.S. patent application Ser. No.14/192,179 entitled “MANAGING SPECULATIVE MEMORY ACCESS REQUESTS IN THEPRESENCE OF TRANSACTIONAL STORAGE ACCESSES,” filed on Feb. 27, 2014, thedisclosure of which is incorporated herein by reference in its entiretyfor all purposes.

BACKGROUND OF THE INVENTION

The present invention relates generally to data processing and, inparticular, to storage accesses to the distributed shared memory systemof a data processing system.

A conventional multiprocessor (MP) computer system, such as a servercomputer system, includes multiple processing units all coupled to asystem interconnect, which typically comprises one or more address, dataand control buses. Coupled to the system interconnect is a systemmemory, which represents the lowest level of volatile memory in themultiprocessor computer system and which generally is accessible forread and write access by all processing units. In order to reduce accesslatency to instructions and data residing in the system memory, eachprocessing unit is typically further supported by a respectivemulti-level cache hierarchy, the lower level(s) of which may be sharedby one or more processor cores.

Cache memories are commonly utilized to temporarily buffer memory blocksthat might be accessed by a processor in order to speed up processing byreducing access latency introduced by having to load needed data andinstructions from system memory. In some MP systems, the cache hierarchyincludes at least two levels. The level one (L1) or upper-level cache isusually a private cache associated with a particular processor core andcannot be accessed by other cores in an MP system. Typically, inresponse to a memory access instruction such as a load or storeinstruction, the processor core first accesses the directory of theupper-level cache. If the requested memory block is not found in theupper-level cache, the processor core then accesses lower-level caches(e.g., level two (L2) or level three (L3) caches) or system memory forthe requested memory block. The lowest level cache (e.g., L3 cache) isoften shared among several processor cores.

In such systems, multiprocessor software concurrently accesses shareddata structures from multiple software threads. When concurrentlyaccessing shared data it is typically necessary to prevent so-called“unconstrained races” or “conflicts”. A conflict occurs between twomemory accesses when they are to the same memory location and at leastone of them is a write and there is no means to ensure the ordering inwhich those accesses occur.

Multiprocessor software typically utilizes lock variables to coordinatethe concurrent reading and modifying of locations in memory in anorderly conflict-free fashion. A lock variable is a location in memorythat is read and then set to a certain value, possibly based on thevalue read, in an atomic fashion. The read-modify-write operation on alock variable is often accomplished utilizing anatomic-read-modify-write (ARMW) instruction or by a sequence ofinstructions that provide the same effect as a single instruction thatatomically reads and modifies the lock variable.

In this manner, a software thread reading an initial “unlocked” valuevia an ARMW instruction is said to have “acquired” the lock and will,until it releases the lock, be the only software thread that holds thelock. The thread holding the lock may safely update the shared memorylocations protected by the lock without conflict with other threadsbecause the other threads cannot obtain the lock until the currentthread releases the lock. When the shared locations have been readand/or modified appropriately, the thread holding the lock releases thelock (e.g., by writing the lock variable to the “unlocked” value) toallow other threads to access the shared locations in storage.

While locking coordinates competing threads' accesses to shared data,locking suffers from a number of well known shortcomings. These include,among others, (1) the possibility of deadlock when a given thread holdsmore than one lock and prevents the forward progress of other threadsand (2) the performance cost of lock acquisition when the lock may nothave been strictly necessary because no conflicting accesses would haveoccurred to the shared data.

To overcome these limitations, the notion of transactional memory can beemployed. In transactional memory, a set of load and/or storeinstructions are treated as a “transaction.” A transaction succeeds whenthe constituent load and store operations can occur atomically without aconflict with another thread. The transaction fails in the presence of aconflict with another thread and can then be re-attempted. If atransaction continues to fail, software may fall back to using lockingto ensure the orderly access of shared data.

To support transactional memory, the underlying hardware tracks thestorage locations involved in the transaction—the transactionfootprint—as the transaction executes for conflicts. If a conflictoccurs in the transaction footprint, the transaction is aborted andpossibly restarted. Use of transactional memory reduces the possibilityof deadlock due to a thread holding multiple locks because, in thetypical case, no locks are held (the transaction simply attempts to makeone or more storage accesses and restarts if a conflict occurs).Further, the processing overhead of acquiring a lock is generallyavoided.

BRIEF SUMMARY

In at least some embodiments, a cache memory of a data processing systemreceives a speculative memory access request including a target addressof data speculatively requested for a processor core. In response toreceipt of the speculative memory access request, transactional memorylogic determines whether or not the target address of the speculativememory access request hits a store footprint of a memory transaction. Inresponse to determining that the target address of the speculativememory access request hits a store footprint of a memory transaction,the transactional memory logic causes the cache memory to rejectservicing the speculative memory access request.

BRIEF DESCRIPTION OF THE SEVERAL VIEWS OF THE DRAWINGS

FIG. 1 is a high-level block diagram of an exemplary data processingsystem in accordance with one embodiment;

FIG. 2 is a more detailed block diagram of an exemplary processing unitin accordance with one embodiment;

FIG. 3 is a detailed block diagram of lower level cache supportingmemory transactions in accordance with one embodiment;

FIG. 4 is an illustrative example of a memory transaction in accordancewith one embodiment;

FIG. 5 is an illustrative example of a memory transaction including asuspended region in accordance with one embodiment;

FIG. 6A depicts execution of an exemplary program illustrating causalityin a multiprocessor data processing system;

FIG. 6B illustrates execution of an exemplary program including memorytransactions to ensure causality;

FIG. 6C depicts execution of an exemplary program including bothtransactional and non-transactional memory accesses;

FIG. 7 illustrates a multiprocessor data processing system including atleast three processor cores that execute the exemplary program of FIG.6C;

FIG. 8 is a high level logical flowchart of an exemplary method by whicha multiprocessor data processing system ensures causality in executionof a program including both transactional and non-transactional memoryaccesses;

FIG. 9 is an illustrative example of a rewind-only memory in accordancewith one embodiment;

FIG. 10 is more detailed view of transactional memory tracking logic inaccordance with one embodiment;

FIG. 11 is a high level logical flowchart of an exemplary method bywhich a rewind-only transaction is processed in accordance with oneembodiment;

FIGS. 12A-12B together form a high level logical flowchart of anexemplary method by which collisions between memory access requests andmemory transactions are detected and resolved in accordance with oneembodiment; and

FIG. 13 is a data flow diagram illustrating a design process.

DETAILED DESCRIPTION

With reference now to the figures, wherein like reference numerals referto like and corresponding parts throughout, and in particular withreference to FIG. 1, there is illustrated a high level block diagramdepicting an exemplary data processing system 100 in accordance with oneembodiment. In the depicted embodiment, data processing system 100 is acache coherent symmetric multiprocessor (SMP) data processing systemincluding multiple processing nodes 102 a, 102 b for processing data andinstructions. Processing nodes 102 are coupled to a system interconnect110 for conveying address, data and control information. Systeminterconnect 110 may be implemented, for example, as a busedinterconnect, a switched interconnect or a hybrid interconnect.

In the depicted embodiment, each processing node 102 is realized as amulti-chip module (MCM) containing four processing units 104 a-104 d,each preferably realized as a respective integrated circuit. Theprocessing units 104 within each processing node 102 are coupled forcommunication to each other and system interconnect 110 by a localinterconnect 114, which, like system interconnect 110, may beimplemented, for example, with one or more buses and/or switches. Systeminterconnect 110 and local interconnects 114 together form a systemfabric.

As described below in greater detail with reference to FIG. 2,processing units 104 each include a memory controller 106 coupled tolocal interconnect 114 to provide an interface to a respective systemmemory 108. Data and instructions residing in system memories 108 cangenerally be accessed, cached and modified by a processor core in anyprocessing unit 104 of any processing node 102 within data processingsystem 100. System memories 108 thus form the lowest level of volatilestorage in the distributed shared memory system of data processingsystem 100. In alternative embodiments, one or more memory controllers106 (and system memories 108) can be coupled to system interconnect 110rather than a local interconnect 114.

Those skilled in the art will appreciate that SMP data processing system100 of FIG. 1 can include many additional non-illustrated components,such as interconnect bridges, non-volatile storage, ports for connectionto networks or attached devices, etc. Because such additional componentsare not necessary for an understanding of the described embodiments,they are not illustrated in FIG. 1 or discussed further herein. Itshould also be understood, however, that the enhancements describedherein are applicable to cache coherent data processing systems ofdiverse architectures and are in no way limited to the generalized dataprocessing system architecture illustrated in FIG. 1.

Multiprocessor data processing system such as data processing system 100of FIG. 1 implement a memory consistency model that specifies the legalpossible executions of a given multiprocessor program with respect tomemory accesses (e.g., among other things, the values that may bereturned by load instructions, the order of writes to memory, thoseinstruction execution dependencies that affect the ordering of memoryaccesses, and the final values for memory locations at the conclusion ofa multiprocessor program). A memory consistency model is specified bytwo major characteristics: ordering of memory access operations andatomicity of store operations.

The ordering of memory operations specifies how memory operations may,if at all, be re-ordered relative to the order of their respective loadand store instructions in the individual threads of execution in themultiprocessor program. Memory consistency models must define orderingof memory access operations in four general cases: (1) ordering of thememory operations for a load instruction to a following loadinstruction, (2) ordering of the memory operations for a loadinstruction to a following store instruction, (3) ordering of the memoryoperations for a store instruction to a following store instruction, and(4) ordering of the memory operations for a store instruction to afollowing load instruction. Strong consistency memory models will, ingeneral, preserve all or at least most of these orderings. Inparticular, many strong consistency memory models enforce the firstthree orderings, but do not enforce store-to-load ordering. Weakconsistency memory models will generally not enforce most or all ofthese orderings.

Atomicity of store operations refers to whether or not a given thread ofexecution can read the value of its own store operation before otherthreads, and furthermore, whether the value written to the distributedshared memory system by the store operation becomes visible to otherthreads in a logically instantaneous fashion or whether the value canbecome visible to other threads at different points in time. A memoryconsistency model is called “multi-copy atomic” if the value written bya store operation of one thread becomes visible to all other threads ina logically instantaneous fashion. In general, strong consistency memorymodels are multi-copy atomic, and weak consistency memory models do notenforce multi-copy atomicity.

In a given multiprocessor program, program semantics often require thatmulti-copy atomicity and/or the various orderings between memory accessoperations are respected. Therefore, in a data processing system 100having a distributed shared memory system that implements a weakconsistency memory model, so called “barrier” (e.g., SYNC) instructionsare typically provided to allow the programmer to specify what memoryaccess operation orderings and atomicity are to be applied duringexecution of the multiprocessor program.

Referring now to FIG. 2, there is depicted a more detailed block diagramof an exemplary processing unit 104 in accordance with one embodiment.In the depicted embodiment, each processing unit 104 is an integratedcircuit including two or more processor cores 200 a, 200 b forprocessing instructions and data. In a preferred embodiment, eachprocessor core 200 is capable of independently executing multiplehardware threads of execution simultaneously. However, in the followingdescription, unless the interaction between threads executing on a sameprocessor core is relevant in a particular context, for simplicity,terms “processor core” and “thread executing on a processor core” areused interchangeably. As depicted, each processor core 200 includes oneor more execution units, such as load-store unit (LSU) 202, forexecuting instructions. The instructions executed by LSU 202 includememory access instructions that request load or store access to a memoryblock in the distributed shared memory system or cause the generation ofa request for load or store access to a memory block in the distributedshared memory system. Memory blocks obtained from the distributed sharedmemory system by load accesses are buffered in one or more registerfiles (RFs) 208, and memory blocks updated by store accesses are writtento the distributed shared memory system from the one or more registerfiles 208.

The operation of each processor core 200 is supported by a multi-levelvolatile memory hierarchy having at its lowest level a shared systemmemory 108 accessed via an integrated memory controller 106, and at itsupper levels, one or more levels of cache memory, which in theillustrative embodiment include a store-through level one (L1) cache 226within and private to each processor core 200, and a respective store-inlevel two (L2) cache 130 for each processor core 200 a, 200 b. In orderto efficiently handle multiple concurrent memory access requests tocacheable addresses, each L2 cache 130 can be implemented with multipleL2 cache slices, each of which handles memory access requests for arespective set of real memory addresses.

Although the illustrated cache hierarchies includes only two levels ofcache, those skilled in the art will appreciate that alternativeembodiments may include additional levels (L3, L4, etc.) of on-chip oroff-chip, private or shared, in-line or lookaside cache, which may befully inclusive, partially inclusive, or non-inclusive of the contentsthe upper levels of cache.

Each processing unit 104 further includes an integrated and distributedfabric controller 216 responsible for controlling the flow of operationson the system fabric comprising local interconnect 114 and systeminterconnect 110 and for implementing the coherency communicationrequired to implement the selected cache coherency protocol. Processingunit 104 further includes an integrated I/O (input/output) controller214 supporting the attachment of one or more I/O devices (not depicted).

In operation, when a hardware thread under execution by a processor core200 includes a memory access instruction requesting a specified memoryaccess operation to be performed, LSU 202 executes the memory accessinstruction to determine the target address (e.g., an effective address)of the memory access request. After translation of the target address toa real address, L1 cache 226 is accessed utilizing the target address.Assuming the indicated memory access cannot be satisfied solely byreference to L1 cache 226, LSU 202 then transmits the memory accessrequest, which includes at least a transaction type (ttype) (e.g., loador store) and the target real address, to its affiliated L2 cache 130for servicing.

It will also be appreciated by those skilled in the art that accesslatency can be improved by prefetching data that is likely to accessedby a processor core 200 into one or more levels of the associated cachehierarchy in advance of need. Accordingly, processing unit 104 caninclude one or more prefetch engines, such as prefetch engine (PFE) 212,that generate prefetch load requests based on historical demand accesspatterns of processor cores 200. Prefetch load requests generated by PFE212 are inherently speculative in that they are generated prior to, andin expectation of, a subsequent demand access request generated throughthe execution of a memory access instruction by LSU 202.

With reference now to FIG. 3, there is illustrated a more detailed blockdiagram of an exemplary embodiment of a lower level cache (e.g., an L2cache 130) that supports memory transactions in accordance with oneembodiment. As shown in FIG. 3, L2 cache 130 includes a cache array 302and a directory 308 of the contents of cache array 302. Although notexplicitly illustrated, cache array 302 preferably is implemented with asingle read port and single write port to reduce the die area requiredto implement cache array 302.

Assuming cache array 302 and directory 308 are set-associative as isconventional, memory locations in system memories 108 are mapped toparticular congruence classes within cache array 302 utilizingpredetermined index bits within the system memory (real) addresses. Theparticular memory blocks stored within the cache lines of cache array302 are recorded in cache directory 308, which contains one directoryentry for each cache line. While not expressly depicted in FIG. 3, itwill be understood by those skilled in the art that each directory entryin cache directory 308 includes various fields, for example, a tag fieldthat identifies the real address of the memory block held in thecorresponding cache line of cache array 302, a state field that indicatethe coherency state of the cache line, an LRU (Least Recently Used)field indicating a replacement order for the cache line with respect toother cache lines in the same congruence class, and inclusivity bitsindicating whether the memory block is held in the associated L1 cache226.

L2 cache 130 includes multiple (e.g., 16) Read-Claim (RC) machines 312for independently and concurrently servicing load (LD) and store (ST)requests received from the affiliated processor core 200. In order toservice remote memory access requests originating from processor cores200 other than the affiliated processor core 200, L2 cache 130 alsoincludes multiple snoop machines 311. Each snoop machine 311 canindependently and concurrently handle a remote memory access request“snooped” from local interconnect 114. As will be appreciated, theservicing of memory access requests by RC machines 312 may require thereplacement or invalidation of memory blocks within cache array 302.Accordingly, L2 cache 130 also includes CO (castout) machines 310 thatmanage the removal and writeback of memory blocks from cache array 302.

L2 cache 130 further includes an arbiter 305 that controls multiplexersM1-M2 to order the processing of local memory access requests and memorytransaction requests (corresponding to the tbegin, tbegin_rot, tend,tabort, and tcheck instructions described further herein) received fromthe affiliated processor core 200 and remote requests snooped on localinterconnect 114. Such requests, including local load and store andmemory transaction requests and remote load and store requests, areforwarded in accordance with the arbitration policy implemented byarbiter 305 to dispatch logic, such as a dispatch pipeline 306, whichprocesses each read/load and store request with respect to directory 308and cache array 302. As described further below, transactional memory(TM) logic 380 processes memory transaction requests and tracks memoryaccess operations within memory transactions to ensure completion of thememory access operations in an atomic manner or to abort the memorytransactions in the presence of conflicts.

L2 cache 130 also includes an RC queue 320 and a CPI (castout pushintervention) queue 318 that respectively buffer data being insertedinto and removed from the cache array 302. RC queue 320 includes anumber of buffer entries that each individually correspond to aparticular one of RC machines 312 such that each RC machine 312 that isdispatched retrieves data from only the designated buffer entry.Similarly, CPI queue 318 includes a number of buffer entries that eachindividually correspond to a particular one of the castout machines 310and snoop machines 311, such that each CO machine 310 and each snooper311 that is dispatched retrieves data from only the respectivedesignated CPI buffer entry.

Each RC machine 312 also has assigned to it a respective one of multipleRC data (RCDAT) buffers 322 for buffering a memory block read from cachearray 302 and/or received from local interconnect 114 via reload bus313. The RCDAT buffer 322 assigned to each RC machine 312 is preferablyconstructed with connections and functionality corresponding to thememory access requests that may be serviced by the associated RC machine312. RCDAT buffers 322 have an associated store data multiplexer M4 thatselects data bytes from among its inputs for buffering in the RCDATbuffer 322 in response unillustrated select signals generated by arbiter305.

In operation, a processor core 200 transmits store requests comprising atransaction type (ttype), target real address and store data to a storequeue (STQ) 304. From STQ 304, the store data are transmitted to storedata multiplexer M4 via data path 324, and the transaction type andtarget address are passed to multiplexer M1. Multiplexer M1 alsoreceives as inputs processor load requests from processor core 200 anddirectory write requests from RC machines 312. In response tounillustrated select signals generated by arbiter 305, multiplexer M1selects one of its input requests to forward to multiplexer M2, whichadditionally receives as an input a remote request received from localinterconnect 114 via remote request path 326. Arbiter 305 scheduleslocal and remote memory access requests for processing and, based uponthe scheduling, generates a sequence of select signals 328. In responseto select signals 328 generated by arbiter 305, multiplexer M2 selectseither the local request received from multiplexer M1 or the remoterequest snooped from local interconnect 114 as the next memory accessrequest to be processed.

The request selected for processing by arbiter 305 is placed bymultiplexer M2 into dispatch pipeline 306. Dispatch pipeline 306preferably is implemented as a fixed duration pipeline in which each ofmultiple possible overlapping requests is processed for a predeterminednumber of clock cycles (e.g., 4 cycles). During the first cycle ofprocessing within dispatch pipeline 306, a directory read is performedutilizing the request address to determine if the request address hitsor misses in directory 308, and if the memory address hits, thecoherency state of the target memory block. The directory information,which includes a hit/miss indication and the coherency state of thememory block, is returned by directory 308 to dispatch pipeline 306 in asubsequent cycle. As will be appreciated, no action is generally takenwithin an L2 cache 130 in response to miss on a remote memory accessrequest; such remote memory requests are accordingly discarded fromdispatch pipeline 306. However, in the event of a hit or miss on a localmemory access request or a hit on a remote memory access request, L2cache 130 will service the memory access request, which for requeststhat cannot be serviced entirely within processing unit 104, may entailcommunication on local interconnect 114 via fabric controller 216.

At a predetermined time during processing of the memory access requestwithin dispatch pipeline 306, arbiter 305 transmits the request addressto cache array 302 via address and control path 330 to initiate a cacheread of the memory block specified by the request address. The memoryblock read from cache array 302 is transmitted via data path 342 toError Correcting Code (ECC) logic 344, which checks the memory block forerrors and, if possible, corrects any detected errors. For processorload requests, the memory block is also transmitted to load datamultiplexer M3 via data path 340 for forwarding to the affiliatedprocessor core 200.

At the last cycle of the processing of a memory access request withindispatch pipeline 306, dispatch pipeline 306 makes a dispatchdetermination based upon a number of criteria, including (1) thepresence of an address collision between the request address and aprevious request address currently being processed by a castout machine310, snoop machine 311 or RC machine 312, (2) the directory information,and (3) availability of an RC machine 312 or snoop machine 311 toprocess the memory access request. If dispatch pipeline 306 makes adispatch determination that the memory access request is to bedispatched, the memory access request is dispatched from dispatchpipeline 306 to an RC machine 312 or a snoop machine 311. If the memoryaccess request fails dispatch, the failure is signaled to the requestor(e.g., local or remote processor core 200) by a retry response. Therequestor may subsequently retry the failed memory access request, ifnecessary.

While an RC machine 312 is processing a local memory access request, theRC machine 312 has a busy status and is not available to service anotherrequest. While an RC machine 312 has a busy status, the RC machine 312may perform a directory write to update the relevant entry of directory308, if necessary. In addition, the RC machine 312 may perform a cachewrite to update the relevant cache line of cache array 302. Directorywrites 408 a, 408 b and cache writes may be scheduled by arbiter 305during any interval in which dispatch pipeline 306 is not alreadyprocessing other requests according to the fixed scheduling of directoryreads and cache reads. When all operations for the given request havebeen completed, the RC machine 312 returns to an unbusy state.

Associated with RC machines 312 is data handling circuitry, differentportions of which are employed during the servicing of various types oflocal memory access requests. For example, for a local load request thathits in directory 308, an uncorrected copy of the target memory block isforwarded from cache array 302 to the affiliated processor core 200 viadata path 340 and load data multiplexer M3 and additionally forwarded toECC logic 344 via data path 342. In the case of an ECC error in thetarget memory block obtained by the local load request, corrected datais forwarded to RCDAT buffer 322 via data path 346 and store datamultiplexer M4 and then from RCDAT 322 to affiliated processor core 200via data path 360 and load data multiplexer M3. For a local storerequest, store data is received within RCDAT buffer 322 from STQ 304 viadata path 324 and store data multiplexer M4, the store is merged withthe memory block read into RCDAT buffer 322 from cache array 302 via ECClogic 344 and store data multiplexer M4, and the merged store data isthen written from RCDAT buffer 322 into cache array 302 via data path362. In response to a local load miss or local store miss, the targetmemory block acquired through issuing a memory access operation on localinterconnect 114 is loaded into cache array 302 via reload bus 313,store data multiplexer M4, RCDAT buffer 322 (with store merge for astore miss) and data path 362.

Referring now to FIG. 4, an illustrative example of a memory transactionis depicted. Those skilled in the art will recognize that the particularsemantics and instructions utilized to implement the various memorytransactions described herein are but some of the numerous possibleimplementations and that the disclosed techniques of implementingtransactional memory are not dependent on the specific instructions andinstruction semantics employed.

Illustrative memory transaction 400 begins at tbegin instruction 402.Tbegin instruction 402 initiates memory transaction 400, causes theprocessor core 200 executing tbegin instruction 402 to take a checkpoint210 of the architected register state of processor core 200, and (e.g.,through a corresponding tbegin request sent to the affiliated L2 cache130) invokes tracking of load and store instructions within thetransaction body 406 to ensure they complete in an atomic fashion orthat memory transaction 400 fails in the presence of a conflict. Memorytransaction 400 additionally includes a branch instruction 404immediately following tbegin instruction 402. When memory transaction400 first executes, the condition code register in processor core 200upon which branch instruction 404 depends is initialized to a value thatcauses the program branch indicated by branch instruction 404 not to betaken and the flow of execution to continue to transaction body 406. Asdiscussed below, in response to failure of memory transaction 400, thecondition code register is set to a different value, and branchinstruction 404 causes execution to branch to a fail handler routine.

In the exemplary embodiment depicted in FIG. 3, TM logic 380 trackstransactional memory access (e.g., load and store) instructions withintransaction body 406 to ensure that they complete in an atomic fashionor that memory transaction 400 fails in the presence of a conflict. Inparticular, TM tracking logic 381 within TM logic 380 includes a numberof entries that indicate which cache lines in cache array 302 areincluded in the transaction footprint (as described below, for example,with reference to FIG. 10). The transaction footprint includes twoportions: the load footprint corresponding to cache lines touched solelyby loads within transaction body 406 (e.g., the cache line at address Ain exemplary memory transaction 400) and the store footprintcorresponding to cache lines touched solely by store instructions or byboth load and store instructions in transaction body 406 (e.g., thecache line at address B in exemplary memory transaction 400).

As further shown in FIG. 3, TM logic 380 further includes transactionalcontrol logic 382, which controls the sequencing of a memory transactionand provides a pass/fail indication 384 and an optional TM killedindication 385 to the associated processor core 200. Pass/failindication 384 indicates to processor core 200 whether or not the memorytransaction successfully committed to the distributed shared memorysystem at the execution of the tend instruction 408 at the end of memorytransaction 400. TM killed indication 385 indicates to processor core200 whether or not a conflict has occurred during the transaction. Inresponse to transactional control logic 382 asserting TM killedindication 385, processor core 200 may, as a performance optimization,optionally abort and restart memory transaction 400 prior to reachingtend instruction 408.

In response to pass/fail indication 384 (or optionally TM killedindication 385) indicating that a conflict has occurred during executionof memory transaction 400, processor core 200 re-establishes itsarchitected register state from the checkpoint 210 taken at theexecution of tbegin instruction 402, invalidates the tentativelymodified cache lines in the store footprint, releases tracking logic381, sets the condition code register such that branch instruction 404will be taken, and transfers control to branch instruction 404. Inaddition, processor core 200 sets a transaction failure cause register(not shown) in processor core 200 to indicate the cause of the memorytransaction's failure. The fail handler routine invoked by branchinstruction 404 may choose to re-attempt memory transaction 400 or fallback to more conventional locking mechanisms, optionally based on thecontent of the transaction failure cause register.

During the execution of a memory transaction, the values stored to thedistributed shared memory system by transaction body 406 (i.e., those inthe store footprint of the memory transaction) are visible only to thethread of the processor core 200 executing the memory transaction.Threads running on the same or other processor cores 200 will not “see”these values until and only if the memory transaction successfullycommits.

For a memory transaction to successfully commit, the load and storeinstructions in transaction body 406 must complete in an atomic fashion(i.e., there must be no conflicts for the cache lines in the memorytransaction's load and store footprints) and the effects of the storeinstructions in transaction body 406 must propagate to all processingunits 104 in data processing system 100 and invalidate any cached copiesof those cache lines held in other processing units 104. If both ofthese conditions hold when tend instruction 408 is executed,transactional control logic 382 indicates to processor core 200 viapass/fail indication 384 that memory transaction 400 passed and commitsall stores performed in transaction body 406 to L2 cache 130, thusmaking them visible to all other threads and processor cores 200 in thesystem simultaneously.

In the following discussion, a load or store instruction will be called“transactional” if that load or store instruction occurs within thetransaction body 406 of a memory transaction 400. Similarly, a load orstore will be called “non-transactional” if it occurs outside atransaction body 406. In one exemplary embodiment, a conflict policy ofdata processing system 100 defines a conflict with another processorcore's memory access to occur for a given memory transaction in any oneof several possible cases. In a first case, a conflict occurs if anon-transactional store from another processor core 200 hits a cacheline within either the given memory transaction's load or storefootprint. In a second case, a conflict occurs if a transactional storefrom another processor core 200 hits a cache line within the givenmemory transaction's load footprint. In a third case, a conflict occursif a non-transactional load hits a cache line within the given memorytransaction's store footprint. In a fourth case, the given memorytransaction has a conflict if one of its transactional loads hits anaddress already extant in the store footprint of another processorcore's memory transaction. In a fifth case, the given memory transactionhas a conflict if one of its transactional stores hits an addressalready extant in the store footprint of another processor core's memorytransaction. The above conflict policy biases in favor of transactionalstores over transactional loads, while allowing transactional andnon-transactional loads to freely intermingle. This exemplary conflictpolicy, a variant of which is described in greater detail below withreference to FIGS. 12A-12B, is but one of several possible embodiments.

With reference now to FIG. 5, there is illustrated a representativememory transaction 500 containing a suspended region. As can be seen bycomparison of FIGS. 4 and 5, memory transaction 500 includes a tbegininstruction 502, branch instruction 504, transaction body 506 and tendinstruction 508, which correspond to tbegin instruction 402, branchinstruction 404, transaction body 406 and tend instruction 408 describedabove. In addition, memory transaction 500 includes a tsuspendinstruction 510 that initiates the start of a suspended region 512. Whena memory transaction is suspended through execution of tsuspendinstruction 510, the load and store footprints currently established forthe enclosing memory transaction containing suspended region 512 remainin place and continue to be tracked by TM tracking logic 381 forconflicts. However, any load or store instructions within suspendedregion 512 are treated as non-transactional loads and stores and followexisting semantics for such loads and stores. In particular, storeswithin suspended region 512 are non-transactional and will commit andbegin propagating to other processors unconditionally. If a store withinsuspended region 512 hits either the load or the store footprint of theenclosing memory transaction, a conflict occurs (which also destroys thetentative transactional version of the cache line in the storefootprint) and is logged by transactional control logic 382. However,this conflict is not acted on until the enclosing memory transaction isresumed upon execution of tresume instruction 514, at which point theprocessor core 200 passes control to branch instruction 504 asdescribed. If a non-transactional load instruction within suspendedregion 512 hits a cache line within the store footprint of the enclosingmemory transaction 500, that load instruction returns the tentativelyupdated value written by a transactional store within the transactionbody 506 unless that value has been overwritten by a non-transactionalstore either by another processor core 200 or by a non-transactionalstore in suspended region 512, in which case the non-transactional loadinstruction returns the current value of the target location.

Use of a suspended region 512 allows the temporary suspension of amemory transaction, which permits store instructions in the suspendedregion 512 to unconditionally update locations in the distributed sharedmemory system while also allowing for the resumption of the memorytransaction at a later time. One possible use for a suspended region 512is to log debug information into a scratchpad region of the distributedshared memory system and then to resume the enclosing memorytransaction. Without a suspended region, the write of the debuginformation would be rolled back any time the enclosing memorytransaction is aborted.

Referring now to FIG. 6A, the execution of an exemplary programillustrating the property of causality in a multiprocessor dataprocessing system is shown. As used herein “causality,” which isdesirable property in multiprocessor programs, is defined as beingpreserved if, during execution of a multiprocessor program, a giventhread of execution cannot read the effects of a computation before thewrites that caused the computation can be read by the given thread.

In the simplified example given in FIG. 6A (as well as those discussedbelow with reference to FIGS. 6B-6C), a multiprocessor program isexecuted by three processor cores 200 of data processing system 100,labeled for ease of reference as processor core 0, processor core 1 andprocessor core 2. In FIG. 6A, processor core 0 executes a storeinstruction 600 that writes a value of 1 to address A in the distributedshared memory system. This update of address A propagates to processorcore 1, and load instruction 610 executed by processor core 1 thereforereturns a value of 1. Even though the memory update made by storeinstruction 600 has propagated to processor core 1, that memory updatemay not yet have propagated to processor core 2. If store instruction614 executes on processor 1 and the associated memory update propagatesto processor 2 before the memory update of store instruction 600propagates to processor 2, causality would be violated because the storeof the value of 1 to address B, which is an effect of the store toaddress A, would be visible to processor core 2 before the memory updateassociated with causal store instruction 600 was visible to processorcore 2.

To ensure causality in a weak consistency memory model, barrierinstruction 612 (e.g., a SYNC) ensures that store instruction 614 doesnot take effect or begin propagating its memory update to otherprocessor cores until load instruction 610 has bound to its value. Inaddition, barrier instruction 612 also ensures that the memory updateassociated with store instruction 600 propagates to processor 2 beforethe memory update associated with store instruction 614. Thus, causalityis preserved because the cause of the computation (i.e., the memoryupdate of store instruction 600) is visible to processor core 2 beforethe result of the computation (i.e., the memory update of store 614). Abarrier instruction 622 is also executed by processor core 2 to ensurethat processor core 2 executes load instructions 620 and 624 and bindstheir values in order, thus guaranteeing that processor core 2 properlyobserves the memory updates made by processor core 0 and processor core1.

With reference now to FIG. 6B, an exemplary embodiment of themultiprocessor program of FIG. 6A rendered in terms of memorytransactions is illustrated. In FIG. 6B, the branch instructions to thememory transaction fail handler are omitted for clarity.

As illustrated, processor core 0 executes a memory transaction 630including a tbegin instruction 632, tend instruction 636, and atransaction body including a store instruction 634 that stores a valueof 1 to address A. Upon the execution of tend instruction 636, memorytransaction 600 successfully commits and makes the update to address Avisible to all the other processor cores simultaneously. In particular,by the time load instruction 642 of the memory transaction 640 executingon processor core 1 can read the value of 1 from address A, loadinstruction 654 of the memory transaction 650 executing on processorcore 2 must also be able to read the value of 1 for address A. Memorytransaction 640 then reads the value of 1 for address A, stores a valueof 1 to address B and successfully commits. Finally, load instruction652 of memory transaction 650 reads a value of 1 for address B, andgiven that memory transaction 640 read a value of 1 for A, loadinstruction 654 must also read a value of 1 for address A.

In order to make the memory updates of store instructions in asuccessful transaction visible to all other processor coressimultaneously, before that memory transaction can commit all the cacheline invalidates necessitated by the memory transaction must havepropagated through the data processing system such that any otherprocessor cores' now stale copies of the updated cache lines have beenremoved (e.g., invalidated) and can no longer be read by the otherprocessor cores. Without this requirement, a processor core could stillread a stale value for an updated memory location after the memorytransaction that updated the memory location committed. A processorcore, therefore, needs to ensure that the memory updates associated withits own transactional stores are fully propagated through the dataprocessing system to invalidate any stale cached copies beforecommitting a successful memory transaction in order to maintain thesemantics of memory transactions. As a consequence of the propagation ofthe memory updates inherent in the semantics of memory transactions,causality is trivially preserved when only memory transactions areutilized to access memory locations in a distributed shared memorysystem. However, when transactional and non-transactional code interacton the same shared variables, causality is not directly preserved byensuring that the memory updates made by a memory transaction arevisible simultaneously to all other processor cores.

Referring now to FIG. 6C, an illustrative multiprocessor program isdepicted that includes a mixture of transactional and non-transactionalaccesses to a distributed shared memory system. In FIG. 6C, the branchinstructions to the memory transaction fail handler are again omittedfor clarity.

In the exemplary multiprocessor program, processor core 0 executes anon-transactional store instruction 660 that unconditionally writes avalue of 1 to address A in the distributed shared memory system. Thisvalue propagates to processor core 1 and is read by transactional loadinstruction 672 within the memory transaction 670 executed by processorcore 1. Processor core 1 then executes a store instruction 674 withinmemory transaction 670 that updates the cache line associated withaddress B and completes invalidating any stale cached copies of thecache line associated with address B (so that no other processor coreholds a copy of the now stale cache line) and successfully commitsmemory transaction 670 upon execution of tend instruction 676. Processorcore 2 then executes load instructions 680 and 684 to read, in order,the cache lines associated with addresses B and A, respectively, basedon the ordering enforced by barrier instruction 682. If transaction 670only ensures that its own memory updates are fully propagated throughthe distributed shared memory system before committing, the memoryupdate of store instruction 660 may or may not have propagated toprocessor core 2. Therefore, in at least some operating scenarios,processor core 2 could read a value of 1 for the cache line associatedwith address B and the, now stale, initial value of 0 for the cache lineassociated with address A, thus violating causality. The same resultwould be obtained if processor core 2 utilized transactional loads toread from addresses A and B, as depicted for processor 2 in FIG. 6B.

To guarantee causality, memory transaction 670 must ensure not only thatits own transactional stores are propagated throughout the entiredistributed shared memory system, but also that any non-transactionalstore that is read by a transactional load within the transaction hasalso propagated throughout the distributed shared memory system. (Memoryupdates of transactional writes that are read by the memory transactionare guaranteed to have propagated throughout the distributed sharedmemory system because those memory updates could not be read bytransaction 670 before they were visible to the entire distributedshared memory system). To ensure that the memory updates ofnon-transactional stores read by memory transaction 670 are alsopropagated throughout the distributed shared memory system, theprocessing of the tend instruction 676 of memory transaction 670 mustnot allow commitment of memory transaction 670 until the memory updateof any non-transactional store read by memory transaction 670 ispropagated throughout the distributed shared memory system.

With reference now to FIG. 7, there is illustrated a partial view ofdata processing system 100 of FIG. 1, which executes the multiprocessorprogram of FIG. 6C. In the view given in FIG. 7, processor cores 200 a,200 b and 200 c respectively correspond to processor cores 0, 1 and 2 ofFIG. 6C. Further, an instance of causality resolution logic 379 isinstantiated for and coupled to each instance of snooper 311, forexample, as a component of the L2 cache 130 affiliated with eachprocessor core 200.

Initially, processor core 200 c holds a cached copy of the initial value(e.g., 0) of memory location A in its L1 cache 226 c. Processor 200 abegins execution of the multiprocessor program of FIG. 6C by executingstore instruction 660. In response to execution of store instruction660, processor core 200 a transmits a store request to its L2 cache 130a, which allocates an RC machine 312 to service the store request. RCmachine 312 broadcasts the store request onto local interconnect 114,and snoop machine 311 c of the L2 cache 130 c affiliated with processorcore 200 c registers the store request, including the processing unitthat sourced the store request (i.e., the processing unit includingprocessor core 200 a). At this point, the memory update of storeinstruction 660 has not propagated to processor core 200 c, but isinstead queued for later processing, advantageously allowing processorcore 200 a to continue executing further instructions before the memoryupdate of store instruction 660 is fully propagated.

Processor core 200 b then executes load instruction 672 and, finding nocopy of the target cache line associated with address A in its L1 cache226 b, transmits a read request to its L2 cache 130 b. In response tothe read request, L2 cache 130 b allocates RC machine 312 b to servicethe read request. In response to a miss of the read request in L2 cache130 b, RC machine 312 b issues a read request onto local interconnect114 to obtain the current value for address A. L2 cache 130 a respondsto the read request and provides the current value of address A toprocessor core 200 b by cache-to-cache intervention. At this point aso-called “causality passing read” has occurred, that is, loadinstruction 672 has read the value of a store instruction that has notfully propagated through the entire distributed shared memory system. Toaccount for this fact and to protect causality, causality resolutionlogic 379 c in L2 cache 130 c notes the successful read interventionbetween the vertical cache hierarchies of processor cores 200 a and 200b for an address that is currently being invalidated by snoop machine311 c. In this manner causality resolution logic 379 c directly tracksthe causal dependency that processor 200 b and its vertical cachehierarchy has on the memory update of store instruction 660 completingits propagation.

Processor 200 b executes store instruction 674, which specifies anupdate of the value of address B to 1. In response to execution of storeinstruction 674, RC machine 312 b issues a store request correspondingto store instruction 674 on local interconnect 114. In absence of anexisting cached copy of the target cache line, memory controller 106supplies the current value of address B from system memory 108 inresponse to the store request, and RC machine 312 b updates L2 cache 130b accordingly. At this point processor core 1 executes tend instruction676 to attempt to successfully commit transaction 670 and places acorresponding TEND request on local interconnect 114 to ensure that allprior memory updates by transactional stores in memory transaction 670have been propagated throughout the distributed shared memory system andthat any memory updates by non-transactional stores read by memorytransaction 670 have similarly propagated throughout the distributedshared memory system. In this case, the memory update of storeinstruction 674 has fully propagated throughout the distributed sharedmemory system because no other caches held a copy of the cache lineassociated with address B. However, had any such copy existed and hadthe memory update not been fully complete, a snoop machine 311 in thosecaches, which noted the initial processor core 200 issuing the store,would be active and would provide a retry response to the snooped TENDrequest from that processor core 200 (forcing the TEND request to bereissued) until the invalidation of the cached copy of the cache linecompletes.

In the case at hand, the TEND request is not from the processor core 200that initiated the store request, and therefore snoop machine 311 c willnot provide a retry response to the TEND request. However, causalityresolution logic 379 c has a causal dependency for processor 200 b andits vertical cache hierarchy and issues on local interconnect 114 aretry response to the TEND request because the TEND request was issuedfrom a processor core 200 that was the recipient of a causality passingread of the same address that snoop machine 311 c is processing. In thismanner, causality resolution logic 379 directly tracks which processorcores 200 have a causality dependency due to reading a memory update ofa non-transactional store that was not fully completed for the processorcore with which causality resolution logic 379 is associated.

It should be noted that, in general, causality resolution logic 379 mustmaintain a list capable of representing all the processors cores 200 inthe data processing system to provide causality in cases in which thecausality dependency chain passes through more than one processor core(e.g., a test where a first processor stores a location, a secondprocessor reads that location and then stores a first flag variable, athird processor loads the first flag variable and writes a second flagin a transaction, and then a final thread reads the second flag and thenthe initial location). In such an implementation, a TEND request issuedfrom any processor core with a causal dependency on the target addressbeing invalidated by the snoop machine 311 associated with the instanceof causality resolution logic 379 is retried. In a large SMP, however,such an embodiment can be prohibitive in cost and many implementationsof causality resolution logic 379 only precisely track causal dependencychains of a certain fixed depth (e.g., two or three processors) and inthe presence of longer dependency chains resort to pessimisticallyretrying all TEND requests until the cache line invalidationsnecessitated by the store instruction have completed processing.

To summarize, causality resolution logic is utilized to detect theoccurrence of causal dependency chains, to a depth determined by theembodiment, on a pending store that has not completed processingthroughout the entire distributed shared memory system. These causaldependencies are utilized to stall the completion of TEND requests fromthose processor cores with a causal dependency on the incomplete(pending) stores. In this manner, the memory transaction cannot complete(and therefore make its own stores visible), until the stores the memorytransaction has read (i.e., those in the causal dependency chain of thememory transaction) have first completed throughout the distributedshared memory system. Only after these stores in the memorytransaction's causal dependency chain (and the transactional stores ofthe memory transaction itself, though this is guaranteed by snooper 311instead of causality resolution logic 379) have completed, may the TENDrequest complete, leading to the memory transaction successfullycommitting if no conflicts have occurred during its execution.

In other embodiments, additional causality resolution logic may berequired to ensure the causality of memory operations. For example, inan implementation that contains a write-through L1 cache shared by amultithreaded processor core followed by a shared L2 store queue, it ispossible for different threads (i.e., logically different processorcores from the point of view of software) to read stored values from theL1 cache before these stores have even propagated to the L2 cache, muchless to the entire distributed shared memory system. In such animplementation, the tend instruction must act as a barrier fortransactional stores in the given thread. This behavior ensures that thetransactional stores are propagated to the system interconnect and thenecessary snoop machines 311 so that the tend instruction can ensure,when trying to complete the memory transaction, that all of the cacheline invalidations required by the memory transaction's stores havefully propagated. In addition, the tend instruction must act as abarrier for non-transactional stores that have been (or may have been)read by transactional loads within the transaction. In the simplest (andmost common embodiment), all non-transactional stores within the sharedstore queue are treated as if they have come from a single thread forpurposes of retrying the TEND request.

In this manner, all non-transactional stores from which any transactionhas (or may have) read that have not been fully propagated are broadcastto snoop machines 311 as necessary before a TEND request for anytransaction from that multithreaded processor core is presented on localinterconnect 114. In such an embodiment, snoop machines 311 treat allstores coming from a given multithreaded processor core in a unifiedmanner and will retry any TEND request, as necessary, from that givenmultithreaded processor core regardless of thread. In this embodiment,causality resolution logic 379 is not involved in monitoring theseintra-core dependencies, but instead is utilized solely to managecausality dependencies between multithreaded processor cores.

The exact placement and details of the necessary causality resolutionlogic will vary with the particulars of given embodiment and will beapparent to those skilled in the art given the teachings herein. Ingeneral, at any point where a load may return the value of a store thathas not fully propagated throughout the entire distributed shared memorysystem, if causality is to be preserved a mechanism must be provided toensure that any store with a causal dependency to a different processorcore is noted and that causal dependency delays the processing of a tendinstruction (or other semantic) ending a memory transaction until suchtime as the stores in the causal dependency chain of the memorytransaction have completed propagating.

Referring now to FIG. 8, there is depicted a high level logicalflowchart of the processing of a tend instruction terminating a memorytransaction in accordance with one embodiment. The process begins atblock 800, for example, in response to initiation of execution of a tendinstruction within the LSU 202 of a processor core 200. The process ofFIG. 8 proceeds from block 800 to block 801, which depicts LSU 202ensuring that all prior suspend mode load instructions and all priortransactional load instructions have their values bound. This checkensures the transactional load instructions are present in the memorytransaction's footprint and that the suspend mode load instructions haveobtained their values. The process proceeds from block 801 to block 802,which depicts ensuring that the cache line invalidations necessitated bytransactional stores within the memory transaction have been fullypropagated throughout the distributed shared memory system. In theembodiment described above, verification of propagation of the cacheline invalidations necessitated by transactional stores is accomplishedby one or more snoop machines 311 providing a retry response to anyapplicable TEND request on local interconnect 114 until the previoustransactional stores have invalidated all cached copies of the memorylocation(s) targeted by the memory updates. The process then proceeds tostep 804, which illustrates ensuring that the cache line invalidationsnecessitated by causally dependent non-transactional stores havecompletely propagated throughout the distributed shared memory system.In the embodiment described above, verification of propagation of thecache line invalidations necessitated by non-transactional stores isaccomplished by one or more instances of causality resolution logic 379providing a retry response to any applicable TEND request on localinterconnect 114 until the previous memory updates of causally dependentnon-transactional stores have invalidated all cached copies of thememory location(s) targeted by the memory updates.

At block 806, transactional control logic 382 determines whether or nota conflict has occurred for the memory transaction. In response totransactional control logic 382 determining that a conflict hasoccurred, the process proceeds to block 808, which depicts transactionalcontrol logic 382 invalidating the tentative store footprint of thememory transaction (e.g., as recorded in L2 cache 130) and indicatingvia pass/fail indication 384 that the memory transaction has failed. Asfurther illustrated at block 808, in response to pass/fail indication384 processor core 200 updates its condition code register and transferscontrol to the fail handling branch instruction within the memorytransaction (block 808). The process then terminates at step 812.

Returning to block 806, in response to transactional control logic 382determining that no conflict has occurred during execution of the memorytransaction, the process proceeds to step 810, which depicts TM controllogic 382 committing the transaction, inter alia, by causing thetransaction footprint to be committed to the distributed shared memorysystem (e.g., by updating one or more coherence states in the directory308 of L2 cache 130 to indicate the transaction footprint is valid andavailable for access by all threads) and indicating to processor core200 via pass/fail indication 384 that the memory transaction passed. Theprocess then terminates at block 812.

Memory transactions, as described above, enable a programmer to enforceexecution of groups of load and/or store instructions by a dataprocessing system in an atomic fashion and to fail and repeat the memorytransactions as necessary to preserve the appearance of atomicity of thestorage accesses of the memory transactions in the presence of conflictswith other storage accesses. While memory transactions provide avaluable and needed capability, there is also a need to be able tospeculatively execute a block of instructions, particularly includingstore instructions, and then to be able to discard the results of thatexecution under software control without regard to the existence ofconflicting accesses. For example, some programming models require thatthe execution of certain code sequences do not cause a fault. To avoidsuch faults, an additional code sequence is typically required tovalidate that the inputs to the code sequence will not produce a faultbefore the sequence is executed. This pre-validation can incursignificant additional overhead. However, with a “rewind only”transaction (ROT) as described herein, the code sequence may bespeculatively executed without the additional overhead of validation andmay then be rewound if a fault occurs.

Discarding or “rewinding” the storage-modifying effects of a storeinstruction has traditionally not been supported in prior processors,and therefore the amount of speculation permitted for a storeinstruction (and for instructions dependent on that store instruction)was severely limited. As described herein, the mechanisms supportingtransactional memory may be adapted, reused and extended to efficientlysupport a discardable speculative execution mechanism for blocks ofinstructions, specifically those including store instructions. Withoutthe enhancements described herein, a full memory transaction would berequired to rewind store instructions, at additional cost as describedbelow.

To support rewinding the storage-modifying effects of storeinstructions, a distinct type of memory transaction referred to hereinas a “rewind only” transaction (ROT) is introduced. Unlike a traditionalmemory transaction, a ROT, by definition, does not require any conflictdetection or atomicity guarantees, but rather only provides a semanticto enforce the discarding of the execution results of a group of one ormore speculatively executed instructions that may include one or morestore instructions. Furthermore, the commitment of a ROT does not dependupon or require the propagation of the invalidations of causallydependent non-transactional stores through the distributed shared memorysystem, as described above with reference to block 804 of FIG. 8.

While conflict detection is not required for the semantic definition ofa ROT, a typical implementation will provide conflict tracking for storeinstructions within the ROT, if only to avoid additional unnecessarycomplexity in the design of the processor core and cache hierarchy atlittle additional benefit. So while conflict tracking is not requiredfor a ROT as a matter of definition (because atomicity is not preservedby a ROT), as a matter of implementation, hardware supporting executionof ROTs will typically provide conflict tracking for the store footprintof a ROT for simplicity.

The utility of retaining store footprint conflict tracking for ROTs canbe seen in the management of conflicts between different threads on amulti-threaded processor core sharing a common write-through L1 cache.In such a multi-threaded processor core, if multiple threads wereconcurrently executing ROTs including store instructions targeting agiven cache line, the L1 cache would have to be able to maintain adifferent image of the given cache line for each thread (i.e., the L1cache would have to be able to hold multiple concurrently active imagesof any given cache line). Furthermore, when each ROT committed, the L1cache would have to be able to merge the updates made to the cache lineby the thread committing the ROT into the remaining concurrent copy orcopies of the cache line—an operation that is exceptionally complex. Ingeneral, it is far more efficient and less costly to employ the existingconflict tracking mechanisms for the store footprint of a ROT as if itwere a non-ROT memory transaction.

Typically, load instructions will significantly outnumber storeinstructions in a memory transaction. For TM control logic 380 of agiven capacity, a significantly larger transaction can therefore beaccommodated as a ROT rather than a non-ROT memory transaction (which,in the absence of ROTs would have to be employed to rewind speculativelyexecuted store instructions). Furthermore, a ROT can successfullycomplete in the presence of false sharing conflicts (i.e., a conflictthat occurs, for example, when a store instruction from another threadwrites within a cache line in the footprint of a memory transaction, butdoes not actually alter the data being manipulated by the memorytransaction). Because conflicts are tracked on a per cache-line basisand not on a per-location basis, such false sharing conflicts cause thefailure of memory transactions that are not strictly required by thedefinition of a memory transaction, but must occur due to thelimitations of the conflict tracking implementation. ROTs, however, aremore resilient in the presence of such false sharing conflicts than thenon-ROT memory transactions that would have to be used in the absence ofsupport for ROTs.

With reference now to FIG. 9, an illustrative example of arepresentative ROT 900 is illustrated. ROT 900 may form, for example, aportion of a multiprocessor program.

ROT 900 begins with a unique instruction, tbegin_rot 902, whichidentifies the beginning of a ROT. Similar to a normal (i.e., non-ROT)memory transaction 400 of FIG. 4, the instruction immediately followingtbegin_rot instruction 902 is a branch instruction 904 that redirectsexecution to a failure handling routine in response to the ROT 900either failing or (as explained below) aborting under software control.Branch instruction 904 is followed by a transaction body 906, which maycontain transactional memory access (e.g., load and/or store) or otherinstructions, and possibly one or more tabort instruction(s) 910. Ifpresent, tabort instruction 910 directs execution of ROT 900 to beaborted and execution results of ROT 900 to be discarded. Although notillustrated in FIG. 9, ROT 900 may further optionally include bypassinstructions that determine if ROT 900 should be aborted (e.g., based ona variable value read from the distributed shared memory system by atransactional load of the ROT or the availability of a system resource)and that, responsive to the determination, either cause tabortinstruction 910 to be executed or cause execution to branch aroundtabort instruction 910 to one or more transactional instructionincluding tend instruction 908, which, when executed, causes ROT 900 tobe committed (and, in particular, makes the storage-modifying effects ofthe store instructions within transaction body 906 non-speculative).

If a tabort instruction 910 within a ROT 900 is executed, the executionresults of ROT 900 are discarded, a condition code register is updatedto indicate that a tabort instruction 910 caused the ROT to fail, andcontrol passes to branch instruction 904, which is taken based on thevalue present in the condition code register. Execution of a tabortinstruction 910 is the primary way in which the speculative executionresults of a ROT 900 are discarded and control is passed to the failhandler via branch instruction 904. Among other reasons, a ROT 900 (ornon-ROT memory transaction) may also fail and pass control to the failhandler via branch instruction 904 (or branch instruction 404) due to acapacity overflow (overflowing the capacity of TM logic 380) or due toexecution of an instruction (e.g., a cache-inhibited load or storeinstruction) that can have untracked side effects and therefore isinherently unable to be re-executed and consequently cannot legallyappear in a ROT or memory transaction (which may have to be executedseveral times to successfully commit).

Referring now to FIG. 10, there is illustrated a more detailed view ofTM tracking logic 381 in accordance with one embodiment. As depicted, TMtracking logic 381 includes a TM directory 1000, which contains a numberof entries 1002 for tracking the cache lines within the load and storefootprints of ROTs and/or non-ROT memory transactions. In the depictedembodiment, each entry 1002 within TM directory 1000 includes threefields: address tag field 1004, load valid (LV) field 1006, and storevalid (SV) field 1008. Address tag field 1004 indicates the real memoryaddress of a cache line that is in the footprint of the ROT or non-ROTmemory transaction. SV field 1006 and LV field 1008 respectivelyindicate whether the cache line is part of the store footprint or loadfootprint of the memory transaction. In at least one embodiment, LVfield 1006 and SV field 1008 are mutually exclusive, meaning that, for agiven entry 1002, one or neither of LV field 1006 and SV field 1008 maybe set concurrently but not both. When both of fields 1006 and 1008 arereset, the entry 1002 is invalid and no cache line is then being trackedby that entry 1002.

For a non-ROT memory transaction, when a transactional load is presentedto TM logic 380 and there is no entry in TM directory 1000 for thetarget cache line of the transactional load, a new entry 1002 isallocated (possibly evicting an existing entry 1002), the address tagfield 1004 of the new entry is updated with the address tag of thetarget cache line, and the LV field 1006 is set. If, on the other hand,an existing entry 1002 is already tracking the target cache line (andtherefore either LV field 1006 or SV field 1008 is already set), noupdate to the existing entry 1002 is made because the target cache lineof the transactional load is already being tracked.

As with a transactional load, if a transactional store of a non-ROTmemory transaction is presented to TM logic 380 and there is no entry inTM directory 1000 for the target cache line of the transactional store,a new entry 1002 is allocated (possibly evicting an existing entry1002), the address tag field 1004 of the new entry is updated with theaddress tag of the target cache line, and the SV field 1008 is set. If,on the other hand, an existing entry 1002 is already tracking the targetcache line and LV field 1006 is set for that entry 1002, then LV field1006 is reset, and SV field 1008 is set to indicate that this cache lineis now part of the store footprint for the memory transaction. If SVfield 1008 is already set for the existing entry 1002, no update toentry 1002 is performed.

In response to a ROT or non-ROT memory transaction committing orfailing, TM tracking logic 381 clears the entries 1002 in TM directory1000.

For a ROT, TM tracking logic 381 updates TM directory 1000 as describedabove for transactional stores of non-ROT memory transactions. However,for loads within the ROT, TM tracking logic 381 does not update TMdirectory 1000 because the load footprint is not tracked for conflictsin a ROT. This behavior can be implemented in at least two ways. In afirst implementation, all non-transactional load and store operationstransmitted from a processor core 200 to its L2 cache 130 are identifiedas either being a non-transactional load or store, a non-ROTtransactional load or store, or as a ROT transactional load or store. Inthis case, TM tracking logic 381 ignores ROT transactional loads forpurposes of updating TM directory 1000. In another implementation, allnon-transactional loads and stores and ROT loads are identified as beingnon-transactional and are accordingly ignored by TM tracking logic 381for purposes of updating TM directory 1000. ROT transactional stores andnon-ROT transactional loads and stores are identified as beingtransactional, and TM tracking logic 381 accordingly updates TMdirectory 1000 as described above for non-ROT transactional loads andstores. In either implementation, TM tracking logic 381 preferably doesnot update TM directory 1000 for ROT transactional loads.

With reference now to FIG. 11, there is illustrated a high level logicalflowchart of a method of processing of a tend instruction terminating anon-ROT memory transaction or a ROT in accordance with one embodiment.For ease of understanding, like reference numerals are utilized todenote steps corresponding to those depicted in FIG. 8.

The process of FIG. 11 begins at block 800, for example, in response toinitiation of execution of a tend instruction within the LSU 202 of aprocessor core 200. LSU 202 then ensures at block 801 that all priorsuspend mode load instructions and all prior non-ROT transactional loadinstructions have their values bound. This check ensures the non-ROTtransactional load instructions are present in the memory transaction'sfootprint and that the suspend mode load instructions have obtainedtheir values. LSU 202 then determines at block 1100 whether or not thememory transaction terminated by the tend is a ROT or non-ROT memorytransaction. In response to a determination that the memory transactionis a non-ROT memory transaction, the process continues to block 802 andsubsequent blocks, which have been described.

Returning to block 1100, in response to a determination that the tendinstruction terminates a ROT, blocks 802, 804 and 806 are bypassed asunnecessary for a ROT, and control passes to block 1102. Block 1102depicts LSU 202 querying TM logic 380 whether a conflict for the ROT'sstore footprint was detected by TM tracking logic 381 (as opposed to aconflict on either the load or store footprint for a non-ROT memorytransaction). In response to TM logic 380 indicating a conflict has beendetected for the store footprint of the ROT, the process proceeds toblock 808, which depicts failing the ROT and invalidating its storefootprint as described above. In response to TM logic 380 indicating atblock 1102 that no conflict for the ROT has been detected, the processproceeds to block 810, which illustrates commitment of the ROT to thedistributed shared memory system as described above. It should again benoted that commitment of the ROT does not require observance ofcausality, as described above with reference to block 804. Followingeither block 808 or block 810, the process concludes at block 812.

Commercially available processors commonly allow out-of-order executionof the instructions within a thread in order to achieve high instructionthroughput. In such processors, numerous instructions within a samethread can be at various stages of execution at any given time, and theprocessor can generally execute the instructions in any order as long asdependencies between instructions (and barrier instructions) areobserved. Commercially available processors also commonly supportspeculative instruction execution, that is, the execution ofinstructions for which it is not certain that the execution results willbe committed to the architected state of the processor. It should beappreciated that out of the various types of memory access instructions,load instructions can be executed speculatively, but store instructions(except for execution within an enclosing memory transaction) generallycannot be executed speculatively because store instructions necessarilymodify the architected state of the processor. The definition of whichload instructions are “speculative” can also vary betweenimplementations. For example, in some implementations, a loadinstruction can be considered speculative if it follows in program ordera branch instruction that has not yet been resolved as taken ornot-taken. In other implementations, a more expansive definition can beemployed. For example, any load instruction can be deemed speculative aslong as it is not the “next to complete” (NTC) instruction, where theNTC instruction is defined as the instruction for which all precedinginstructions in program order have been architecturally committed. Asemployed herein, however, the term “speculative” is defined to excludethose instructions whose execution results may or may not be committedto the architected state of the processor only by virtue of inclusion ofthose instructions within a memory transaction. (Transactionalinstructions may, however, still be speculative, as made clear below.)

In the absence of memory transactions, speculative load instructions canbe discarded and re-executed one or many times without any negativearchitectural side effects and with a relatively small performanceimpact. However, if speculative load instructions are executed in thepresence of memory transactions, speculative load requests generated bythe execution of speculative load instructions of a first thread can hitthe footprint of a memory transaction of a second thread tracked by TMtracking logic 381 and cause such memory transactions to fail. Suchcollisions can occur both in cases in which the first and second threadsare executed on different processor cores 200 and in cases in which thefirst and second threads are simultaneously executed on a commonprocessor core 200 and share a common vertical cache hierarchy.

The present disclosure recognizes that any memory transaction failurescaused by speculative load instructions that are ultimately discardedare unnecessary from an architectural perspective and amplify thenegative performance impact of unsuccessful speculative execution. Dataprefetch requests, which are inherently speculative, can also have theundesirable effect of unnecessarily causing memory transactions to fail.To reduce unnecessary failures of memory transactions, TM logic, such asTM logic 380, preferably responds to speculative memory access requestsof one thread (e.g., speculative load requests and data prefetchrequests) that hit the store footprint of a memory transaction ofanother thread with a response and/or special data packet that causesthe speculative memory access request to fail and be re-executed, ifnecessary, rather than the memory transaction failing.

Referring now to FIGS. 12A-12B, there is depicted a high level logicalflowchart of an exemplary method by which collisions between memoryaccess requests and memory transactions are detected and resolved inaccordance with one embodiment. The illustrated method can be performed,for example, by TM logic 380 of FIG. 3.

The process begins at block 1200, for example, in response to receipt byTM logic 380 of a memory access request, which can be received from theassociated processor core 200 or received from another processor core200 via local interconnect 114. In response to receipt of the memoryaccess request, TM logic 380 determines at block 1202 whether or not thememory access request is speculative. As noted above, store requests arenon-speculative, but all data prefetch requests and at least some loadrequests, whether transactional or non-transactional, are speculative.In at least one embodiment, load requests defined to be speculative inthe relevant embodiment of data processing system 100 are identified assuch by a ttype or label associated with the load requests by theissuing processor core 200.

In response to a determination at block 1202 that the memory accessrequest received by TM logic 380 is not a speculative request (i.e., isa store request or a non-speculative load request), the process passesthrough page connector A to block 1230 of FIG. 12B, which is describedbelow. If, on the other hand, TM logic 380 determines at block 1202 thatthe memory access request is a speculative request (i.e., a load requestor data prefetch request), TM logic 380 further determines at block 1204whether or not the speculative memory access request is transactional,that is, generated by execution of a load instruction within a memorytransaction. In at least one embodiment, TM logic 380 determines whetheror not the memory access request is transactional, for example, by attype or label associated with the memory access request. In response toa determination at block 1204 that the memory access request istransactional and therefore a transactional load request (and not a dataprefetch request), the process proceeds to block 1212, which isdescribed below. In response to determination at block 1204 that thespeculative memory access request is not transactional and is thereforea non-transactional load request or data prefetch request, the processpasses to block 1206.

As indicated at block 1206, if TM logic 380 is not presently configured(e.g., in hardware or by software) to filter speculativenon-transactional load and data prefetch requests, TM logic 380 handlesthe speculative non-transactional load or data prefetch request asdepicted at block 1208 and following blocks. If, on the other hand, TMlogic 380 is presently configured to filter speculativenon-transactional load and data prefetch requests, TM logic 380 handlesthe speculative non-transactional load or data prefetch request asdepicted at block 1214 and following blocks.

Referring now to block 1208, TM logic 380 determines whether or not thetarget address of the speculative non-transactional load or dataprefetch request hits the store footprint of a memory transaction ofanother thread that is being tracked by TM tracking logic 381. If not,TM logic 380 allows the speculative non-transactional load or dataprefetch request to be serviced in accordance with the coherenceprotocol implemented by data processing system 100 (block 1216).Following block 1216, the process of FIG. 12A ends at block 1220. If,however, TM logic 380 determines at block 1208 that the target addressof the speculative non-transactional load or data prefetch request hitsthe store footprint of a memory transaction of another thread that isbeing tracked by TM tracking logic 381, TM logic 380 fails the memorytransaction as shown at block 1210 by returning to the associatedprocessor core 200 a fail indication 384 at the conclusion of the memorytransaction, and optionally, an earlier TM killed indication 385.Thereafter, the process passes to blocks 1216 and 1220, which have beendescribed.

Referring now to block 1212, if TM logic 380 is not presently configured(e.g., in hardware or by software) to filter speculative transactionalmemory access requests (i.e., speculative transactional load requests),TM logic 380 handles the speculative transactional load request asdepicted at block 1208 and following blocks, which have been described.If, on the other hand, TM logic 380 is presently configured to filterspeculative transactional load requests, TM logic 380 handles thespeculative non-transactional load request as depicted at block 1214 andfollowing blocks.

At block 1214, TM logic 380 determines whether or not the target addressof the speculative transactional load request hits the store footprintof a memory transaction of another thread that is being tracked by TMtracking logic 381. If not, TM logic 380 allows the speculativetransactional load request to be serviced in accordance with thecoherence protocol implemented by data processing system 100 (block1216). Thereafter, the process of FIG. 12A ends at block 1220. If,however, TM logic 380 determines at block 1214 that the target addressof the speculative transactional load request hits the store footprintof a memory transaction of another thread that is being tracked by TMtracking logic 381, TM logic 380 causes L2 cache 230 to not service thespeculative transactional memory access request and to provide arejection indication to the requesting processor core 200 as shown atblock 1218. In various embodiments, a rejection indication provided tothe associated processor core 200 can include, for example, a retrysignal and/or an indicator provided with the data. A rejectionindication provided in response to a speculative transactional loadrequest snooped on local interconnect 114 can include, for example, acoherence response message (e.g., a partial response representing theindividual coherence response of L2 cache 230 to the snooped loadrequest) and/or an indicator provided with the data. Following block1218, the process of FIG. 12A ends at block 1220.

It should be noted that in the process of FIG. 12A no check is madewhether the target address of the speculative load or data prefetchrequest hits the load footprint of a memory transaction because anaddress collision between a speculative load or data prefetch requestand the load footprint has no effect on the successful completion of thememory transaction.

Referring now to FIG. 12B, if the memory access request received by TMlogic 380 is non-speculative (i.e., a non-speculative load request or astore request), TM logic 380 determines at block 1230 whether or not thenon-speculative load or store request is transactional. If not, theprocess proceeds to block 1232, which illustrates TM logic 380determining (e.g., by reference to the ttype) whether or not thenon-speculative non-transactional memory access request is a loadrequest. In response to a determination at block 1232 that thenon-speculative, non-transactional memory access request is a loadrequest, TM logic 380 determines whether or not the target address ofthe load request hits the store footprint of the memory transaction ofanother thread (block 1234). If not, TM logic 380 allows thenon-speculative, non-transactional load request to be serviced by L2cache 230 in accordance with the coherence protocol implemented by dataprocessing system 100 (block 1238). Thereafter, the method of FIG. 12Bends at block 1280. If, however, TM logic 380 determines at block 1234that the target address of the non-speculative, non-transactional loadrequest hits the store footprint of the memory transaction of anotherthread, TM logic 380 fails the memory transaction as shown at block 1236by returning a fail indication 384 to the associated processor core 200at the conclusion of the memory transaction, and optionally, an earlierTM killed indication 385. Thereafter, the process passes to blocks 1238and 1280, which have been described.

Referring again to block 1232, in response to a determination that thenon-speculative non-transactional memory access request is a storerequest, TM logic 380 determines whether or not the target address ofthe store request hits the load or store footprint of the memorytransaction of another thread (block 1240). If not, TM logic 380 allowsthe processing of any existing transaction to continue as indicated bythe process passing to block 1280 and terminating. If, however, TM logic380 determines at block 1240 that the target address of the storerequest hits the load or store footprint of the memory transaction ofanother thread, TM logic 380 fails the existing memory transaction asshown at block 1242 by returning to the associated processor core 200 afail indication 384 at the conclusion of the memory transaction, andoptionally, an earlier TM killed indication 385. Thereafter, the processof FIG. 12B ends at block 1280.

Referring again to block 1230, in response to TM logic 380 determiningthat the non-speculative memory access request is transactional, theprocess proceeds to block 1250, which illustrates TM logic 380determining (e.g., by reference to the ttype) whether or not thenon-speculative transactional memory access request is a transactionalload request. In response to a determination at block 1250 that thenon-speculative transactional memory access request is a transactionalload request, TM logic 380 determines whether or not the target addressof the load request hits the store footprint of the memory transactionof another thread (block 1270). If not, TM logic 380 allows thenon-speculative transactional load request to be serviced in accordancewith the coherence protocol implemented by data processing system 100(block 1274). Thereafter, the method of FIG. 12B ends at block 1280. IfTM logic 380 determines at block 1270 that the target address of theload request hits the store footprint of the memory transaction ofanother thread, TM logic 380 also allows the non-speculativetransactional load request to be serviced in accordance with thecoherence protocol implemented by data processing system 100 (block1272). However, TM logic 380 also fails the existing memory transactionas shown at block 1260 by returning to the associated processor core 200a fail indication 384 at the conclusion of the memory transaction, andoptionally, an earlier TM killed indication 385. Thereafter, the processof FIG. 12B ends at block 1280.

Referring again to block 1250, if the non-speculative transactionalmemory access request is a transactional store request, TM logic 380further determines whether or not the target address of thetransactional store request hits the load or store footprint of thememory transaction of another thread, as shown atblocks 1252 and 1254,respectively. In response to a determination that the target address ofthe transactional store request does not hit the load or store footprintof an existing memory transaction of another thread, TM logic 380 allowsthe existing transaction, if any, to proceed as indicated by the processpassing from blocks 1252-1254 to block 1280 and terminating. If,however, TM logic 380 determines at blocks 1252 and 1254 that the targetaddress of the store request hits the load or store footprint of thememory transaction of another thread, TM logic 380 fails either a firstmemory transaction including the transactional store instruction thatgenerated the transactional store request (i.e., the “requestingtransaction”) or a second memory transaction to which the conflictingtransaction footprint belongs (i.e., the “existing transaction”).According to a preferred embodiment of the conflict resolution policywhich favors stores over loads, if TM logic 380 determines at block 1254that the target address of the transactional store request of a first(or “requesting”) transaction hits the store footprint of a second (or“existing”) memory transaction, TM logic 380 fails the first (or“requesting”) transaction as shown at block 1256. On the other hand, ifTM logic 380 determines at block 1252 that the target address of thetransactional store request of a first (or “requesting”) transactionhits the load footprint of a second (or “existing”) memory transaction,TM logic 380 preferably instead fails the second (or “existing”)transaction, as shown at block 1260. As noted above, TM logic 380 failsa transaction at block 1256 or bock 1260 by returning to the associatedprocessor core 200 a fail indication 384 at the conclusion of thetransaction, and optionally, an earlier TM killed indication 385.Following either block 1256 or block 1260, the process of FIG. 12B endsat block 1280.

With reference now to FIG. 13, there is depicted a block diagram of anexemplary design flow 1300 used for example, in semiconductor IC logicdesign, simulation, test, layout, and manufacture. Design flow 1300includes processes, machines and/or mechanisms for processing designstructures or devices to generate logically or otherwise functionallyequivalent representations of the design structures and/or devicesdescribed above and shown in FIGS. 1-3, 7, and 10. The design structuresprocessed and/or generated by design flow 1300 may be encoded onmachine-readable transmission or storage media to include data and/orinstructions that when executed or otherwise processed on a dataprocessing system generate a logically, structurally, mechanically, orotherwise functionally equivalent representation of hardware components,circuits, devices, or systems. Machines include, but are not limited to,any machine used in an IC design process, such as designing,manufacturing, or simulating a circuit, component, device, or system.For example, machines may include: lithography machines, machines and/orequipment for generating masks (e.g. e-beam writers), computers orequipment for simulating design structures, any apparatus used in themanufacturing or test process, or any machines for programmingfunctionally equivalent representations of the design structures intoany medium (e.g. a machine for programming a programmable gate array).

Design flow 1300 may vary depending on the type of representation beingdesigned. For example, a design flow 1300 for building an applicationspecific IC (ASIC) may differ from a design flow 1300 for designing astandard component or from a design flow 1300 for instantiating thedesign into a programmable array, for example a programmable gate array(PGA) or a field programmable gate array (FPGA) offered by Altera® Inc.or Xilinx® Inc.

FIG. 13 illustrates multiple such design structures including an inputdesign structure 1020 that is preferably processed by a design process1310. Design structure 1320 may be a logical simulation design structuregenerated and processed by design process 1310 to produce a logicallyequivalent functional representation of a hardware device. Designstructure 1320 may also or alternatively comprise data and/or programinstructions that when processed by design process 1310, generate afunctional representation of the physical structure of a hardwaredevice. Whether representing functional and/or structural designfeatures, design structure 1320 may be generated using electroniccomputer-aided design (ECAD) such as implemented by a coredeveloper/designer. When encoded on a machine-readable datatransmission, gate array, or storage medium, design structure 1320 maybe accessed and processed by one or more hardware and/or softwaremodules within design process 1310 to simulate or otherwise functionallyrepresent an electronic component, circuit, electronic or logic module,apparatus, device, or system such as those shown in FIGS. 1-3, 7, and10. As such, design structure 1320 may comprise files or other datastructures including human and/or machine-readable source code, compiledstructures, and computer-executable code structures that when processedby a design or simulation data processing system, functionally simulateor otherwise represent circuits or other levels of hardware logicdesign. Such data structures may include hardware-description language(HDL) design entities or other data structures conforming to and/orcompatible with lower-level HDL design languages such as Verilog andVHDL, and/or higher level design languages such as C or C++.

Design process 1310 preferably employs and incorporates hardware and/orsoftware modules for synthesizing, translating, or otherwise processinga design/simulation functional equivalent of the components, circuits,devices, or logic structures shown in FIGS. 1-3, 7, and 10 to generate anetlist 1380 which may contain design structures such as designstructure 1320. Netlist 1380 may comprise, for example, compiled orotherwise processed data structures representing a list of wires,discrete components, logic gates, control circuits, I/O devices, models,etc. that describes the connections to other elements and circuits in anintegrated circuit design. Netlist 1380 may be synthesized using aniterative process in which netlist 1380 is resynthesized one or moretimes depending on design specifications and parameters for the device.As with other design structure types described herein, netlist 1380 maybe recorded on a machine-readable storage medium or programmed into aprogrammable gate array. The medium may be a non-volatile storage mediumsuch as a magnetic or optical disk drive, a programmable gate array, acompact flash, or other flash memory. Additionally, or in thealternative, the medium may be a system or cache memory, or bufferspace.

Design process 1310 may include hardware and software modules forprocessing a variety of input data structure types including netlist1380. Such data structure types may reside, for example, within libraryelements 1330 and include a set of commonly used elements, circuits, anddevices, including models, layouts, and symbolic representations, for agiven manufacturing technology (e.g., different technology nodes, 32 nm,45 nm, 90 nm, etc.). The data structure types may further include designspecifications 1340, characterization data 1350, verification data 1360,design rules 1370, and test data files 1385 which may include input testpatterns, output test results, and other testing information. Designprocess 1310 may further include, for example, standard mechanicaldesign processes such as stress analysis, thermal analysis, mechanicalevent simulation, process simulation for operations such as casting,molding, and die press forming, etc. One of ordinary skill in the art ofmechanical design can appreciate the extent of possible mechanicaldesign tools and applications used in design process 1310 withoutdeviating from the scope and spirit of the invention. Design process1310 may also include modules for performing standard circuit designprocesses such as timing analysis, verification, design rule checking,place and route operations, etc.

Design process 1310 employs and incorporates logic and physical designtools such as HDL compilers and simulation model build tools to processdesign structure 1320 together with some or all of the depictedsupporting data structures along with any additional mechanical designor data (if applicable), to generate a second design structure 1390.Design structure 1390 resides on a storage medium or programmable gatearray in a data format used for the exchange of data of mechanicaldevices and structures (e.g., information stored in a IGES, DXF,Parasolid XT, JT, DRG, or any other suitable format for storing orrendering such mechanical design structures). Similar to designstructure 1320, design structure 1390 preferably comprises one or morefiles, data structures, or other computer-encoded data or instructionsthat reside on transmission or data storage media and that whenprocessed by an ECAD system generate a logically or otherwisefunctionally equivalent form of one or more of the embodiments of theinvention shown in FIGS. 1-3, 7, and 10. In one embodiment, designstructure 1390 may comprise a compiled, executable HDL simulation modelthat functionally simulates the devices shown in FIGS. 1-3, 7, and 10.

Design structure 1390 may also employ a data format used for theexchange of layout data of integrated circuits and/or symbolic dataformat (e.g., information stored in a GDSII (GDS2), GL1, OASIS, mapfiles, or any other suitable format for storing such design datastructures). Design structure 1390 may comprise information such as, forexample, symbolic data, map files, test data files, design contentfiles, manufacturing data, layout parameters, wires, levels of metal,vias, shapes, data for routing through the manufacturing line, and anyother data required by a manufacturer or other designer/developer toproduce a device or structure as described above and shown in FIGS. 1-3,7, and 10. Design structure 1390 may then proceed to a stage 1395 where,for example, design structure 1390: proceeds to tape-out, is released tomanufacturing, is released to a mask house, is sent to another designhouse, is sent back to the customer, etc.

As has been described, in at least one embodiment, a cache memory of adata processing system receives a speculative memory access requestincluding a target address of data speculatively requested for aprocessor core. In response to receipt of the speculative memory accessrequest, transactional memory logic determines whether or not the targetaddress of the speculative memory access request hits a store footprintof a memory transaction. In response to determining that the targetaddress of the speculative memory access request hits a store footprintof a memory transaction, the transactional memory logic causes the cachememory to reject servicing the speculative memory access request.

While various embodiments have been particularly shown and described, itwill be understood by those skilled in the art that various changes inform and detail may be made therein without departing from the spiritand scope of the appended claims and these alternate implementations allfall within the scope of the appended claims. For example, althoughaspects have been described with respect to a computer system executingprogram code that directs the functions of the present invention, itshould be understood that present invention may alternatively beimplemented as a program product including a computer-readable storagedevice storing program code that can be processed by a data processingsystem. The computer-readable storage device can include volatile ornon-volatile memory, an optical or magnetic disk, or the like, butexcludes signal media.

As an example, the program product may include data and/or instructionsthat when executed or otherwise processed on a data processing systemgenerate a logically, structurally, or otherwise functionally equivalentrepresentation (including a simulation model) of hardware components,circuits, devices, or systems disclosed herein. Such data and/orinstructions may include hardware-description language (HDL) designentities or other data structures conforming to and/or compatible withlower-level HDL design languages such as Verilog and VHDL, and/or higherlevel design languages such as C or C++. Furthermore, the data and/orinstructions may also employ a data format used for the exchange oflayout data of integrated circuits and/or symbolic data format (e.g.information stored in a GDSII (GDS2), GL1, OASIS, map files, or anyother suitable format for storing such design data structures).

What is claimed is:
 1. A method of data processing in a data processingsystem, the method comprising: at a cache memory of a data processingsystem, receiving a speculative memory access request, the speculativememory access request including a target address of data speculativelyrequested for a processor core; in response to receipt of thespeculative memory access request, transactional memory logicdetermining whether the target address of the speculative memory accessrequest hits a store footprint of a memory transaction; and in responseto determining that the target address of the speculative memory accessrequest hits a store footprint of a memory transaction, thetransactional memory logic causing the cache memory to reject servicingthe speculative memory access request.
 2. The method of claim 1, whereinthe speculative memory access request is a load request.
 3. The methodof claim 1, wherein the speculative memory access request comprises adata prefetch request.
 4. The method of claim 1, wherein: thedetermining includes determining whether the speculative memory accessrequest is a transactional speculative memory access request or anon-transactional speculative memory access request; causing the cachememory to reject servicing the speculative memory access requestcomprises causing the cache memory to reject servicing the speculativememory access request in response to determining the memory accessrequest is a transactional speculative memory access request; and themethod further comprises failing the memory transaction in response todetermining the speculative memory access request is a non-transactionalspeculative memory access request.
 5. The method of claim 1, wherein:the determining includes determining whether the speculative memoryaccess request is a transactional speculative memory access request or anon-transactional speculative memory access request; causing the cachememory to reject servicing the speculative memory access requestcomprises causing the cache memory to reject servicing the speculativememory access request in response to determining the memory accessrequest is a non-transactional speculative memory access request; andthe method further comprises failing the memory transaction in responseto determining the speculative memory access request is a transactionalspeculative memory access request.
 6. The method of claim 1, wherein:the memory transaction is second memory transaction; the target addressis a second target address; the method further comprises: in response toreceiving at the cache memory a non-speculative transactional loadrequest of a first memory transaction, the non-speculative transactionalload request including a first target address that hits the storefootprint of the second memory transaction, the transactional memorylogic failing the first memory transaction.